On the No-Counterexample Interpretation

1999 ◽  
Vol 64 (4) ◽  
pp. 1491-1511 ◽  
Author(s):  
Ulrich Kohlenbach
Keyword(s):  
Normal Form ◽  
Formal System ◽  
Modus Ponens ◽  
Peano Arithmetic ◽  
Order Type ◽  
Cut Elimination ◽  
Substitution Method ◽  
Functional Interpretation ◽  
First Order ◽  
Definition Of

AbstractIn [15], [16] G. Kreisel introduced the no-counterexample interpretation (n.c.i.) of Peano arithmetic. In particular he proved, using a complicated ε-substitution method (due to W. Ackermann), that for every theorem A (A prenex) of first-order Peano arithmetic PA one can find ordinal recursive functionals of order type < ε0 which realize the Herbrand normal form AH of A.Subsequently more perspicuous proofs of this fact via functional interpretation (combined with normalization) and cut-elimination were found. These proofs however do not carry out the no-counterexample interpretation as a local proof interpretation and don't respect the modus ponens on the level of the nocounterexample interpretation of formulas A and A → B. Closely related to this phenomenon is the fact that both proofs do not establish the condition (δ) and—at least not constructively—(γ) which are part of the definition of an ‘interpretation of a formal system’ as formulated in [15].

scholarly journals On the No-Counterexample Interpretation

1997 ◽  
Vol 4 (42) ◽  
Author(s):  
Ulrich Kohlenbach
Keyword(s):  
Normal Forms ◽  
Formal System ◽  
Modus Ponens ◽  
Order Type ◽  
Cut Elimination ◽  
Gamma Delta ◽  
Functional Interpretation ◽  
First Order ◽  
Epsilon Substitution ◽  
Definition Of

In [15],[16] Kreisel introduced the no-counterexample interpretation (n.c.i.) of Peano<br />arithmetic. In particular he proved, using a complicated epsilon-substitution method (due to<br />W. Ackermann), that for every theorem A (A prenex) of first-order Peano arithmetic (PA) one can find ordinal recursive functionals Phi_A of order type < epsilon_0 which realize the<br />Herbrand normal form A^H of A.<br /> Subsequently more perspicuous proofs of this fact via functional interpretation (combined<br />with normalization) and cut-elimination were found. These proofs however do not carry out the n.c.i. as a local proof interpretation and don't respect the modus<br />ponens on the level of the n.c.i. of formulas A and A -> B. Closely related to this phenomenon is the fact that both proofs do not establish the condition (delta) and - at<br />least not constructively - (gamma) which are part of the definition of an `interpretation of a<br />formal system' as formulated in [15].<br />In this paper we determine the complexity of the n.c.i. of the modus ponens rule for<br />(i) PA-provable sentences,<br />(ii) for arbitrary sentences A;B in L(PA) uniformly in functionals satisfying the n.c.i. of (prenex normal forms of) A and A -> B; and<br />(iii) for arbitrary A;B in L(PA) pointwise in given alpha(< epsilon_0)-recursive functionals satisfying the n.c.i. of A and A -> B. <br /> This yields in particular perspicuous proofs of new uniform versions of the conditions ( gamma), (delta). <br /> Finally we discuss a variant of the concept of an interpretation presented in [17] and<br />show that it is incomparable with the concept studied in [15],[16]. In particular we show<br />that the n.c.i. of PA_n by alpha(<omega_n(omega))-recursive functionals (n >= 1) is an interpretation in the sense of [17] but not in the sense of [15] since it violates the condition (delta).

scholarly journals Extracting Herbrand Disjunctions by Functional Interpretation

2003 ◽  
Vol 10 (32) ◽  
Author(s):  
Philipp Gerhardy ◽  
Ulrich Kohlenbach
Keyword(s):  
Normal Form ◽  
Polynomial Complexity ◽  
Predicate Logic ◽  
Extraction Technique ◽  
Cut Elimination ◽  
Functional Interpretation ◽  
First Order ◽  
Term Extraction ◽  
Higher Types ◽  
First Order Predicate Logic

Carrying out a suggestion by Kreisel, we adapt Gödel's functional interpretation to ordinary first-order predicate logic(PL) and thus devise an algorithm to extract Herbrand terms from PL-proofs. The extraction is carried out in an extension of PL to higher types. The algorithm consists of two main steps: first we extract a functional realizer, next we compute the beta-normal-form of the realizer from which the Herbrand terms can be read off. Even though the extraction is carried out in the extended language, the terms are ordinary PL-terms. In contrast to approaches to Herbrand's theorem based on cut elimination or epsilon-elimination this extraction technique is, except for the normalization step, of low polynomial complexity, fully modular and furthermore allows an analysis of the structure of the Herbrand terms, in the spirit of Kreisel, already prior to the normalization step. It is expected that the implementation of functional interpretation in Schwichtenberg's MINLOG system can be adapted to yield an efficient Herbrand-term extraction tool.

A proof of the cut-elimination theorem in simple type theory

1973 ◽  
Vol 38 (2) ◽  
pp. 215-226
Author(s):  
Satoko Titani
Keyword(s):  
Boolean Algebra ◽  
Type Theory ◽  
Arbitrary Number ◽  
Formal System ◽  
Cut Elimination ◽  
Simple Type ◽  
Inductive Definition ◽  
Simple Type Theory ◽  
Maximal Ideals ◽  
Definition Of

In [4], I introduced a quasi-Boolean algebra, and showed that in a formal system of simple type theory, from which the cut rule is omitted, wffs form a quasi-Boolean algebra, and that the cut-elimination theorem can be formulated in algebraic language. In this paper we use the result of [4] to prove the cut-elimination theorem in simple type theory. The theorem was proved by M. Takahashi [2] in 1967 by using the concept of Schütte's semivaluation. We use maximal ideals of a quasi-Boolean algebra instead of semivaluations.The logical system we are concerned with is a modification of Schütte's formal system of simple type theory in [1] into Gentzen style.Inductive definition of types.0 and 1 are types.If τ1, …, τn are types, then (τ1, …, τn) is a type.Basic symbols.a1τ, a2τ, … for free variables of type τ.x1τ, x2τ, … for bound variables of type τ.An arbitrary number of constants of certain types.An arbitrary number of function symbols with certain argument places.

On the definition of ‘formal deduction’

1956 ◽  
Vol 21 (2) ◽  
pp. 129-136 ◽  
Author(s):  
Richard Montague ◽  
Leon Henkin
Keyword(s):  
Functional Calculus ◽  
Modus Ponens ◽  
Formal Proof ◽  
Axiom Schema ◽  
First Order ◽  
Individual Variable ◽  
Individual Variables ◽  
Definition Of ◽  
Formal Rules

The following remarks apply to many functional calculi, each of which can be variously axiomatized, but for clarity of exposition we shall confine our attention to one particular system Σ. This system is to have the usual primitive symbols and formation rules of the pure first-order functional calculus, and the following formal axiom schemata and formal rules of inference.Axiom schema 1. Any tautologous wff (well-formed formula).Axiom schema 2. (a) A ⊃ B, where A is any wff, a and b are any individual variables, and B arises from A by replacing all free occurrences of a by free occurrences of b.Axiom schema 3. (a)(A ⊃ B)⊃(A⊃ (a)B). where A and B are any wffs, and a is any individual variable not free in A.Rule of Modus Ponens: applies to wffs A and A ⊃ B, and yields B.Rule of Generalization: applies to a wff A and yields (a)A, where a is any individual variable.A formal proof in Σ is a finite column of wffs each of whose lines is a formal axiom or arises from two preceding lines by the Rule of Modus Ponens or arises from a single preceding line by the Rule of Generalization. A formal theorem of Σ is a wff which occurs as the last line of some formal proof.

scholarly journals An application of graphical enumeration to PA

2003 ◽  
Vol 68 (1) ◽  
pp. 5-16
Author(s):  
Andreas Weiermann
Keyword(s):  
Normal Form ◽  
Asymptotic Analysis ◽  
Natural Number ◽  
Proof Theory ◽  
Peano Arithmetic ◽  
First Order ◽  
Primitive Recursive Arithmetic ◽  
Independence Result ◽  
Primitive Recursive

AbstractFor α less than ε0 let Nα be the number of occurrences of ω in the Cantor normal form of α. Further let ∣n∣ denote the binary length of a natural number n, let ∣n∣h denote the h-times iterated binary length of n and let inv(n) be the least h such that ∣n∣h ≤ 2. We show that for any natural number h first order Peano arithmetic, PA, does not prove the following sentence: For all K there exists an M which bounds the lengths n of all strictly descending sequences 〈α0, …, αn〉 of ordinals less than ε0 which satisfy the condition that the Norm Nαi of the i-th term αi is bounded by K + ∣i∣ · ∣i∣i.As a supplement to this (refined Friedman style) independence result we further show that e.g., primitive recursive arithmetic, PRA, proves that for all K there is an M which bounds the length n of any strictly descending sequence 〈α0, …, αn〉 of ordinals less than ε0 which satisfies the condition that the Norm Nαi of the i-th term αi is bounded by K + ∣i∣ · inv(i). The proofs are based on results from proof theory and techniques from asymptotic analysis of Polya-style enumerations.Using results from Otter and from Matoušek and Loebl we obtain similar characterizations for finite bad sequences of finite trees in terms of Otter's tree constant 2.9557652856.…

Gödel’s Proof Based on ω-Consistency

1992 ◽  
Author(s):  
Raymond M. Smullyan
Keyword(s):  
Predicate Symbol ◽  
Order Theory ◽  
Modus Ponens ◽  
Axiom System ◽  
Peano Arithmetic ◽  
Order Logic ◽  
First Order Logic ◽  
Underlying Assumption ◽  
First Order ◽  
First Order Theory

The proof that we have just given of the incompleteness of Peano Arithmetic was based on the underlying assumption that Peano Arithmetic is correct—i.e., that every sentence provable in P.A. is a true sentence. Gödel’s original incompleteness proof involved a much weaker assumption—that of ω-consistency to which we now turn. We consider an arbitrary axiom system S whose formulas are those of Peano Arithmetic, whose axioms include all those of Groups I and II (or alternatively, any set of axioms for first-order logic with identity such that all logically valid formulas are provable from them), and whose inference rules are modus ponens and generalization. (It is also possible to axiomatize first-order logic in such a way that modus ponens is the only inference rule—cf. Quine [1940].) In place of the axioms of Groups III and IV, however, we can take a completely arbitrary set of axioms. Such a system S is an example of what is termed a first-order theory, and we will consider several such theories other than Peano Arithmetic. (For the more general notion of a first-order theory, the key difference is that we do not necessarily start with + and × as the undefined function symbols, nor do we necessarily take ≤ as the undefined predicate symbol. Arbitrary function symbols and predicate symbols can be taken, however, as the undefined function and predicate symbols—cf. Tarski [1953] for details. However, the only theories (or “systems”, as we will call them) that we will have occasion to consider are those whose formulas are those of P.A.) S is called simply consistent (or just “consistent” for short) if no sentence is both provable and refutable in S.

The substitution method

1965 ◽  
Vol 30 (2) ◽  
pp. 175-192 ◽  
Author(s):  
W. W. Tait
Keyword(s):  
Free Variable ◽  
Order System ◽  
Conservative Extension ◽  
Substitution Method ◽  
Recursive Definition ◽  
First Order ◽  
Variable Elimination ◽  
Primitive Recursion ◽  
Definition Of ◽  
Primitive Recursive

This paper deals with Hilbert's substitution method for eliminating bound variables from first order proofs. With a first order system S framed in the ε-calculus [2] the problem is to associate a system S' without bound variables and an effective procedure for transforming derivations in S into derivations in S′. The transform of a formula A derived in S is to be an “ε-substitution instance” of A, i.e. it is obtained by replacing terms εxB(x) in A by terms of S′. In general the choice of these terms will depend on the particular derivation of A, and not on A alone. Cf. [4]. The present formulation sharpens Hilbert's original statement of the problem, i.e. that the transform of A should be finitistically verifiable, by making explicit the methods of verification used, namely those formalized in S′; on the other hand, it generalizes Hilbert's formulation since S′ need not be restricted to finitist systems.The bound variable elimination procedure can always be taken to be primitive recursive in (the Gödel number of) the derivation of A. Constructions which transcend primitive recursion can simply be built into S′.In this paper we show that if S′ is taken to be a second order system with constants for functionals, then the existence of suitable ε-substitution instances can be expressed by the solvability of certain functional equations in S′. We deal with two cases here. If S is number theory without induction, i.e. essentially predicate calculus with identity, then we can solve the equations in question by taking for S′ the free variable part S* of S with an added rule of definition of functionals by cases (recursive definition on finite ordinals), which is a conservative extension of S*.

Two interpolation theorems for a predicate calculus

1971 ◽  
Vol 36 (2) ◽  
pp. 262-270
Author(s):  
Shoji Maehara ◽  
Gaisi Takeuti
Keyword(s):  
Normal Form ◽  
Completeness Theorem ◽  
Interpolation Theorem ◽  
Second Order ◽  
Predicate Calculus ◽  
Cut Elimination ◽  
First Order ◽  
Image Position ◽  
Order Predicate Calculus

A second order formula is called Π1 if, in its prenex normal form, all second order quantifiers are universal. A sequent F1, … Fm → G1 …, Gn is called Π1 if a formulais Π1If we consider only Π1 sequents, then we can easily generalize the completeness theorem for the cut-free first order predicate calculus to a cut-free Π1 predicate calculus.In this paper, we shall prove two interpolation theorems on the Π1 sequent, and show that Chang's theorem in [2] is a corollary of our theorem. This further supports our belief that any form of the interpolation theorem is a corollary of a cut-elimination theorem. We shall also show how to generalize our results for an infinitary language. Our method is proof-theoretic and an extension of a method introduced in Maehara [5]. The latter has been used frequently to prove the several forms of the interpolation theorem.

On the nonexistence of certain normal forms in the logic of provability

1982 ◽  
Vol 47 (3) ◽  
pp. 638-640
Author(s):  
George Boolos
Keyword(s):  
Normal Form ◽  
Normal Forms ◽  
Modus Ponens ◽  
Propositional Modal Logic ◽  
Formal Systems ◽  
Normal Form Theorem ◽  
Modal Sentence ◽  
The One ◽  
Definition Of ◽  
Functional Combination

G is the system of propositional modal logic whose axioms are all tautologies and all sentences □(A → B) → (□A → □B) and □(□A → A) → □A and whose rules are modus ponens and necessitation. For the connections between G and provability in formal systems, see [3] and [2].A letterless sentence is a modal sentence that contains no sentence letters at all. In [1] the author showed (in effect) the existence of simple normal forms in G for letterless sentences: every letterless sentence is equivalent in G to a truth-functional compound of sentences ⋄r ⊤ (⋄0B = B; ⋄r+1B = ⋄ ⋄r+1 B. ⊤ is the 0-ary connective that is always evaluated as true; we assume that the language contains ⊤ as a primitive.)A family of sets {Hn}n∈ω of modal sentences that contain no sentence letters other than some fixed one (say p) was introduced by Solovay in [3]: H0 = the set of sentences equivalent in G to one of p, −p, ⊤, and ⊥; Hn+x = the set of sentences equivalent in G to a truth-functional combination of sentences ⋄rB, where r ∈ ω and B ∈ Hn. (This definition of the Hn's differs inessentially from the one given in [3].) Since r may = 0, Hn ⊆ Hn+1 and thus if m ≤ n, Hm ⊆ Hn; every modal sentence containing no letter other than p is in some Hn. It follows from the normal form theorem for letterless sentences that every letterless sentence is in H1.In [3] Solovay stated without proof a result about the sets Hn: for every n, Hn ⊊ Hn+1; equivalently, for every n, there is a sentence containing no letter other than p not in Hn. Thus the existence of a normal form for letterless sentences is a very special feature of these sentences, for Solovay's result shows that normal forms like those obtainable for letterless sentences are not to be found in general. He takes this result to be one reason for regarding the Lindenbaum algebra in G of sentences containing no letter other than p as much more complicated than that of letterless sentences.

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